Concurrency Control [R&G] Chapter 17 CS4320 1 Conflict - - PowerPoint PPT Presentation

CS4320 1

Concurrency Control

[R&G] Chapter 17

CS4320 2

Conflict Serializable Schedules

Two schedules are conflict equivalent if:

Involve the same actions of the same transactions Every pair of conflicting actions is ordered the

same way

Schedule S is conflict serializable if S is

conflict equivalent to some serial schedule

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Example

A schedule that is not conflict serializable: The cycle in the graph reveals the problem.

The output of T1 depends on T2, and vice- versa.

T1: R(A), W(A), R(B), W(B) T2: R(A), W(A), R(B), W(B) T1 T2 A B Dependency graph

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Dependency Graph

Dependency graph: One node per Xact; edge

from Ti to Tj if Tj reads/writes an object last written by Ti.

Theorem: Schedule is conflict serializable if

and only if its dependency graph is acyclic

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Review: Strict 2PL

Strict Two-phase Locking (Strict 2PL) Protocol:

Each Xact must obtain a S (shared) lock on object

before reading, and an X (exclusive) lock on object before writing.

All locks held by a transaction are released when

the transaction completes

• If an Xact holds an X lock on an object, no other

Xact can get a lock (S or X) on that object.

Strict 2PL allows only schedules whose

precedence graph is acyclic

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Two-Phase Locking (2PL)

Two-Phase Locking Protocol

Each Xact must obtain a S (shared) lock on object

before reading, and an X (exclusive) lock on object before writing.

A transaction can not request additional locks

• nce it releases any locks.
• If an Xact holds an X lock on an object, no other

Xact can get a lock (S or X) on that object.

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View Serializability

Schedules S1 and S2 are view equivalent if:

If Ti reads initial value of A in S1, then Ti also reads

initial value of A in S2

If Ti reads value of A written by Tj in S1, then Ti also

reads value of A written by Tj in S2

If Ti writes final value of A in S1, then Ti also writes

final value of A in S2 T1: R(A) W(A) T2: W(A) T3: W(A) T1: R(A),W(A) T2: W(A) T3: W(A)

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Lock Management

Lock and unlock requests are handled by the lock

manager

Lock table entry:

Number of transactions currently holding a lock Type of lock held (shared or exclusive) Pointer to queue of lock requests

Locking and unlocking have to be atomic operations Lock upgrade: transaction that holds a shared lock

can be upgraded to hold an exclusive lock

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Deadlocks

Deadlock: Cycle of transactions waiting for

locks to be released by each other.

Two ways of dealing with deadlocks:

Deadlock detection Deadlock prevention

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Deadlock Detection

Assumption: if a lock request cannot be

satisfied, the transaction is blocked and must wait until the resource becomes available.

Create a waits-for graph:

Nodes are transactions There is an edge from Ti to Tj if Ti is waiting for Tj

to release a lock

Periodically check for cycles in the waits-for

graph

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Deadlock Detection (Continued)

Example: T1: S(A), R(A), S(B) T2: X(B),W(B) X(C) T3: S(C), R(C) X(A) T4: X(B) T1 T2 T4 T3 T1 T2 T3 T3

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Deadlock Prevention

Assign priorities based on timestamps.

Assume Ti wants a lock that Tj holds. Two policies are possible:

Wait-Die: It Ti has higher priority, Ti waits for Tj;

• therwise Ti aborts

Wound-wait: If Ti has higher priority, Tj aborts;

• therwise Ti waits

If a transaction re-starts, make sure it has its

• riginal timestamp

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Multiple-Granularity Locks

Hard to decide what granularity to lock

(tuples vs. pages vs. tables).

Shouldn’t have to decide! Data “containers” are nested:

Tuples Tables Pages Database contains

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Solution: New Lock Modes, Protocol

Allow Xacts to lock at each level, but with a

special protocol using new “intention” locks:

Before locking an item, Xact

must set “intention locks”

• n all its ancestors.

For unlock, go from specific

to general (i.e., bottom-up).

SIX mode: Like S & IX at

the same time.

• IS

IX

• IS

IX √ √ √ √ √ √ S X √ √ S X √ √ √ √ √ √ √ √

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Multiple Granularity Lock Protocol

Each Xact starts from the root of the hierarchy. To get S or IS lock on a node, must hold IS or IX

• n parent node.

What if Xact holds SIX on parent? S on parent?

To get X or IX or SIX on a node, must hold IX or

SIX on parent node.

Must release locks in bottom-up order.

Protocol is correct in that it is equivalent to directly setting locks at the leaf levels of the hierarchy.

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Examples

T1 scans R, and updates a few tuples:

T1 gets an SIX lock on R, then repeatedly gets an S lock on tuples of R, and occasionally upgrades to X on the tuples.

T2 uses an index to read only part of R:

T2 gets an IS lock on R, and repeatedly gets an S lock on tuples of R.

T3 reads all of R:

T3 gets an S lock on R. OR, T3 could behave like T2; can use lock escalation to decide which.

• IS

IX

• IS

IX √ √ √ √ √ √ S X √ √ S X √ √ √ √ √ √ √ √

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Dynamic Databases

If we relax the assumption that the DB is a

fixed collection of objects, even Strict 2PL will not assure serializability:

T1 locks all pages containing sailor records with rating = 1, and finds oldest sailor (say, age = 71). Next, T2 inserts a new sailor; rating = 1, age = 96. T2 also deletes oldest sailor with rating = 2 (and, say, age = 80), and commits. T1 now locks all pages containing sailor records with rating = 2, and finds oldest (say, age = 63).

No consistent DB state where T1 is “correct”!

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The Problem

T1 implicitly assumes that it has locked the

set of all sailor records with rating = 1.

Assumption only holds if no sailor records are added while T1 is executing! Need some mechanism to enforce this

• assumption. (Index locking and predicate

locking.)

Example shows that conflict serializability

guarantees serializability only if the set of

• bjects is fixed!

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Index Locking

If there is a dense index on the rating field

using Alternative (2), T1 should lock the index page containing the data entries with rating = 1.

If there are no records with rating = 1, T1 must lock the index page where such a data entry would be, if it existed!

If there is no suitable index, T1 must lock all

pages, and lock the file/table to prevent new pages from being added, to ensure that no new records with rating = 1 are added.

r=1 Data Index

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Predicate Locking

Grant lock on all records that satisfy some

logical predicate, e.g. age > 2*salary.

Index locking is a special case of predicate

locking for which an index supports efficient implementation of the predicate lock.

What is the predicate in the sailor example?

In general, predicate locking has a lot of

locking overhead.

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Locking in B+ Trees

How can we efficiently lock a particular leaf

node?

Btw, don’t confuse this with multiple granularity locking!

One solution: Ignore the tree structure, just lock

pages while traversing the tree, following 2PL.

This has terrible performance!

Root node (and many higher level nodes) become bottlenecks because every tree access begins at the root.

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Two Useful Observations

Higher levels of the tree only direct searches

for leaf pages.

For inserts, a node on a path from root to

modified leaf must be locked (in X mode, of course), only if a split can propagate up to it from the modified leaf. (Similar point holds w.r.t. deletes.)

We can exploit these observations to design

efficient locking protocols that guarantee serializability even though they violate 2PL.

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A Simple Tree Locking Algorithm

Search: Start at root and go down;

repeatedly, S lock child then unlock parent.

Insert/Delete: Start at root and go down,

• btaining X locks as needed. Once child is

locked, check if it is safe:

If child is safe, release all locks on ancestors.

Safe node: Node such that changes will not

propagate up beyond this node.

Inserts: Node is not full. Deletes: Node is not half-empty.

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Example

ROOT

A B C D E F G H I

20 35 20* 38 44 22* 23* 24* 35* 36* 38* 41* 44* Do: 1) Search 38* 2) Delete 38* 3) Insert 45* 4) Insert 25* 23

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A Better Tree Locking Algorithm (See Bayer-Schkolnick paper)

Search: As before. Insert/Delete:

Set locks as if for search, get to leaf, and set X lock on leaf. If leaf is not safe, release all locks, and restart Xact using previous Insert/Delete protocol.

Gambles that only leaf node will be modified;

if not, S locks set on the first pass to leaf are

• wasteful. In practice, better than previous alg.

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Example

ROOT

A B C D E F G H I

20 35 20* 38 44 22* 23* 24* 35* 36* 38* 41* 44* Do: 1) Delete 38* 2) Insert 25* 4) Insert 45* 5) Insert 45*, then 46* 23

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Even Better Algorithm

Search: As before. Insert/Delete:

Use original Insert/Delete protocol, but set IX locks instead of X locks at all nodes. Once leaf is locked, convert all IX locks to X locks top-down: i.e., starting from node nearest to root. (Top-down reduces chances

• f deadlock.)

(Contrast use of IX locks here with their use in multiple-granularity locking.)

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Hybrid Algorithm

The likelihood that we really need an X lock

decreases as we move up the tree.

Hybrid approach:

Set S locks Set SIX locks Set X locks

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Optimistic CC (Kung-Robinson)

Locking is a conservative approach in which

conflicts are prevented. Disadvantages: Lock management overhead. Deadlock detection/resolution. Lock contention for heavily used objects.

If conflicts are rare, we might be able to gain

concurrency by not locking, and instead checking for conflicts before Xacts commit.

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Kung-Robinson Model

Xacts have three phases:

READ: Xacts read from the database, but make changes to private copies of objects. VALIDATE: Check for conflicts. WRITE: Make local copies of changes public.

ROOT

• ld

new modified

• bjects

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Validation

Test conditions that are sufficient to ensure

that no conflict occurred.

Each Xact is assigned a numeric id.

Just use a timestamp.

Xact ids assigned at end of READ phase, just

before validation begins. (Why then?)

ReadSet(Ti): Set of objects read by Xact Ti. WriteSet(Ti): Set of objects modified by Ti.

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Test 1

For all i and j such that Ti < Tj, check that Ti

completes before Tj begins.

Ti Tj

R V W R V W

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Test 2

For all i and j such that Ti < Tj, check that:

Ti completes before Tj begins its Write phase + WriteSet(Ti) ReadSet(Tj) is empty.

Ti Tj

R V W R V W

Does Tj read dirty data? Does Ti overwrite Tj’s writes?

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Test 3

For all i and j such that Ti < Tj, check that:

Ti completes Read phase before Tj does + WriteSet(Ti) ReadSet(Tj) is empty + WriteSet(Ti) WriteSet(Tj) is empty.

Ti Tj

R V W R V W

Does Tj read dirty data? Does Ti overwrite Tj’s writes?

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Comments on Validation

Assignment of Xact id, validation, and the

Write phase are inside a critical section!

I.e., Nothing else goes on concurrently. If Write phase is long, major drawback.

Optimization for Read-only Xacts:

Don’t need critical section (because there is no Write phase).

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Overheads in Optimistic CC

Must record read/write activity in ReadSet and

WriteSet per Xact.

Must create and destroy these sets as needed.

Must check for conflicts during validation, and

must make validated writes ``global’’.

Critical section can reduce concurrency. Scheme for making writes global can reduce clustering

• f objects.

Optimistic CC restarts Xacts that fail validation.

Work done so far is wasted; requires clean-up.

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``Optimistic’’ 2PL

If desired, we can do the following:

Set S locks as usual. Make changes to private copies of objects. Obtain all X locks at end of Xact, make writes global, then release all locks.

In contrast to Optimistic CC as in Kung-

Robinson, this scheme results in Xacts being blocked, waiting for locks.

However, no validation phase, no restarts (modulo deadlocks).

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Timestamp CC

Idea: Give each object a read-timestamp

(RTS) and a write-timestamp (WTS), give each Xact a timestamp (TS) when it begins: If action ai of Xact Ti conflicts with action aj

• f Xact Tj, and TS(Ti) < TS(Tj), then ai must
• ccur before aj. Otherwise, restart

violating Xact.

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When Xact T wants to read Object O

If TS(T) < WTS(O), this violates timestamp

• rder of T w.r.t. writer of O.

So, abort T and restart it with a new, larger TS. (If restarted with same TS, T will fail again! Contrast use of timestamps in 2PL for ddlk prevention.)

If TS(T) > WTS(O):

Allow T to read O. Reset RTS(O) to max(RTS(O), TS(T))

Change to RTS(O) on reads must be written to

disk! This and restarts represent overheads.

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When Xact T wants to Write Object O

If TS(T) < RTS(O), this violates timestamp order

• f T w.r.t. writer of O; abort and restart T.

If TS(T) < WTS(O), violates timestamp order of

T w.r.t. writer of O.

Thomas Write Rule: We can safely ignore such

• utdated writes; need not restart T! (T’s write is

effectively followed by another write, with no intervening reads.) Allows some serializable but non conflict serializable schedules:

Else, allow T to write O.

T1 T2 R(A) W(A) Commit W(A) Commit

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Timestamp CC and Recoverability

Timestamp CC can be modified

to allow only recoverable schedules: Buffer all writes until writer commits (but update WTS(O) when the write is allowed.) Block readers T (where TS(T) > WTS(O)) until writer of O commits.

Similar to writers holding X locks until commit,

but still not quite 2PL.

T1 T2 W(A) R(A) W(B) Commit

Unfortunately, unrecoverable

schedules are allowed:

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Multiversion Timestamp CC

Idea: Let writers make a “new” copy while

readers use an appropriate “old” copy:

O O’ O’’

MAIN SEGMENT (Current versions of DB objects) VERSION POOL (Older versions that may be useful for some active readers.)

Readers are always allowed to proceed.

– But may be blocked until writer commits.

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Multiversion CC (Contd.)

Each version of an object has its writer’s TS as

its WTS, and the TS of the Xact that most recently read this version as its RTS.

Versions are chained backward; we can

discard versions that are “too old to be of interest”.

Each Xact is classified as Reader or Writer.

Writer may write some object; Reader never will. Xact declares whether it is a Reader when it begins.

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Reader Xact

For each object to be read:

Finds newest version with WTS < TS(T). (Starts with current version in the main segment and chains backward through earlier versions.)

Assuming that some version of every object

exists from the beginning of time, Reader Xacts are never restarted.

However, might block until writer of the appropriate version commits.

T

• ld new

WTS timeline

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Writer Xact

To read an object, follows reader protocol. To write an object:

Finds newest version V s.t. WTS < TS(T). If RTS(V) < TS(T), T makes a copy CV of V, with a pointer to V, with WTS(CV) = TS(T), RTS(CV) = TS(T). (Write is buffered until T commits; other Xacts can see TS values but can’t read version CV.) Else, reject write. T

• ld new

WTS

CV

V

RTS(V)

CS4320 46

Transaction Support in SQL-92

Each transaction has an access mode, a

diagnostics size, and an isolation level.

No No No Serializable Maybe No No Repeatable Reads Maybe Maybe No Read Committed Maybe Maybe Maybe Read Uncommitted Phantom Problem Unrepeatable Read Dirty Read Isolation Level

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Summary

There are several lock-based concurrency

control schemes (Strict 2PL, 2PL). Conflicts between transactions can be detected in the dependency graph

The lock manager keeps track of the locks

• issued. Deadlocks can either be prevented or

detected.

Naïve locking strategies may have the

phantom problem

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Summary (Contd.)

Index locking is common, and affects

performance significantly.

Needed when accessing records via index. Needed for locking logical sets of records (index locking/predicate locking).

Tree-structured indexes:

Straightforward use of 2PL very inefficient. Bayer-Schkolnick illustrates potential for improvement.

In practice, better techniques now known; do

record-level, rather than page-level locking.

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Summary (Contd.)

Multiple granularity locking reduces the overhead

involved in setting locks for nested collections of objects (e.g., a file of pages); should not be confused with tree index locking!

Optimistic CC aims to minimize CC overheads in an

``optimistic’’ environment where reads are common and writes are rare.

Optimistic CC has its own overheads however; most

real systems use locking.

SQL-92 provides different isolation levels that control

the degree of concurrency

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Summary (Contd.)

Timestamp CC is another alternative to 2PL; allows

some serializable schedules that 2PL does not (although converse is also true).

Ensuring recoverability with Timestamp CC requires

ability to block Xacts, which is similar to locking.

Multiversion Timestamp CC is a variant which ensures

that read-only Xacts are never restarted; they can always read a suitable older version. Additional

• verhead of version maintenance.